Method and apparatus for increased performance of a parked data bus in the non-parked direction

ABSTRACT

A distributed system structure for a large-way, symmetric multiprocessor system using a bus-based cache-coherence protocol is provided. The distributed system structure contains an address switch, multiple memory subsystems, and multiple master devices, either processors, I/O agents, or coherent memory adapters, organized into a set of nodes supported by a node controller. The node controller receives transactions from a master device, communicates with a master device as another master device or as a slave device, and queues transactions received from a master device. Since the achievement of coherency is distributed in time and space, the node controller helps to maintain cache coherency. In addition, a bus arbiter in the node controller parks a data bus towards a memory subsystem. The node controller does not use data buffer reservations. The data bus grant line to the memory controller is overloaded to use it as a back-pressure, get-off-the-bus signal as well as a normal data bus grant line. The fairness of the bus is thereby increased by creating a mechanism for getting a “parked” device off the data bus without the use of another dedicated signal between physical components. To ensure that the node controller may stream data to the memory subsystem, the bus is not reparked towards the memory subsystem until a configurable number of cycles after the data bus has been granted to the node controller.

CROSS-REFERENCE TO RELATED APPLICATIONS

The present invention is related to the following applications entitled“METHOD AND APPARATUS FOR PROVIDING GLOBAL COHERENCE IN A LARGE-WAY,HIGH PERFORMANCE SMP SYSTEM”, U.S. application Ser. No. 09/350,032,filed on Jul. 8, 1999; “METHOD AND APPARATUS FOR ACHIEVING CORRECT ORDERAMONG BUS MEMORY TRANSACTIONS IN A PHYSICALLY DISTRIBUTED SMP SYSTEM”,U.S. application Ser. No. 09/350,030, filed on Jul. 8, 1999; “METHOD ANDAPPARATUS USING A DISTRIBUTED SYSTEM STRUCTURE TO SUPPORT BUS-BASEDCACHE-COHERENCE PROTOCOLS FOR SYMMETRIC MULTIPROCESSORS”, U.S.application Ser. No. 09/350,031, filed on Jul. 8, 1999; “METHOD ANDSYSTEM FOR RESOLUTION OF TRANSACTION COLLISIONS TO ACHIEVE GLOBALCOHERENCE IN A DISTRIBUTED SYMMETRIC MULTIPROCESSOR SYSTEM”, U.S.application Ser. No. 09/392,833, filed on Sep. 9, 1999; “METHOD ANDSYSTEM FOR IMPLEMENTING REMSTAT PROTOCOL UNDER INCLUSION ANDNON-INCLUSION OF L1 DATA IN L2 CACHE TO PREVENT READ-READ DEADLOCK”,U.S. application Ser. No. 09/404,400, filed on Sep. 23, 1999; and“METHOD AND SYSTEM FOR CONTROLLING DATA TRANSFERS WITH PHYSICALSEPARATION OF DATA FUNCTIONALITY FROM ADDRESS AND CONTROL FUNCTIONALITYIN A DISTRIBUTED MULTI-BUS MULTIPROCESSOR SYSTEM”, U.S. application Ser.No. 09/404,280, filed on Sep. 23, 1999; “METHOD AND APPARATUS TODISTRIBUTE INTERRUPTS TO MULTIPLE INTERRUPT HANDLERS IN A DISTRIBUTEDSYMMETRIC MULTIPROCESSOR SYSTEM”, U.S. application Ser. No. 09/436,201,filed on Nov. 8, 1999; “METHOD AND APPARATUS TO ELIMINATE FAILED SNOOPSOF TRANSACTIONS CAUSED BY BUS TIMING CONFLICTS IN A DISTRIBUTEDSYMMETRIC MULTIPROCESSOR SYSTEM”, U.S. application Ser. No. 09/436,203,filed on Nov. 8, 1999; “METHOD AND APPARATUS FOR TRANSACTION PACING TOREDUCE DESTRUCTIVE INTERFERENCE BETWEEN SUCCESSIVE TRANSACTIONS IN ADISTRIBUTED SYMMETRIC MULTIPROCESSOR SYSTEM”, U.S. application Ser. No.09/436,204, filed on Nov. 8, 1999; “METHOD AND APPARATUS FOR FAIR DATABUS PARKING PROTOCOL WITHOUT DATA BUFFER RESERVATIONS AT THE RECEIVER”,U.S. application Ser. No. 09/436,202, filed on Nov. 8, 1999; “METHOD ANDAPPARATUS FOR AVOIDING DATA BUS GRANT STARVATION IN A NON-FAIR,PRIORITIZED ARBITER FOR A SPLIT BUS SYSTEM WITH INDEPENDENT ADDRESS ANDDATA BUS GRANTS”, U.S. application Ser. No. 09/436,200, filed on Nov. 8,1999; “METHOD AND APPARATUS FOR SYNCHRONIZING MULTIPLE BUS ARBITERS ONSEPARATE CHIPS TO GIVE SIMULTANEOUS GRANTS FOR THE PURPOSE OF BREAKINGLIVELOCKS”, U.S. application Ser. No. 09/436,192, filed on Nov. 8, 1999;“METHOD AND APPARATUS FOR TRANSACTION TAG ASSIGNMENT AND MAINTENANCE INA DISTRIBUTED SYMMETRIC MULTIPROCESSOR SYSTEM”, U.S. application Ser.No. 09/436,205, filed on Nov. 8, 1999; “METHOD AND SYSTEM FOR DATA BUSLATENCY REDUCTION USING TRANSFER SIZE PREDICTION FOR SPLIT BUS DESIGNS”,U.S. application Ser. No. 09/434,764, filed on Nov. 4, 1999; all ofwhich are assigned to the same assignee.

BACKGROUND OF THE INVENTION

1. Technical Field

The present invention relates generally to an improved data processingsystem and, in particular, to a method and system for improving datathroughput within a data processing system. Specifically, the presentinvention relates to a method and system for improving performance ofinput/output processing and bus access regulation.

2. Description of Related Art

Traditionally, symmetric multiprocessors are designed around a commonsystem bus on which all processors and other devices such as memory andI/O are connected by merely making physical contacts to the wirescarrying bus signals. This common bus is the pathway for transferringcommands and data between devices and also for achieving coherence amongthe system's cache and memory. A single-common-bus design remains apopular choice for multiprocessor connectivity because of the simplicityof system organization.

This organization also simplifies the task of achieving coherence amongthe system's caches. A command issued by a device gets broadcast to allother system devices simultaneously and in the same clock cycle that thecommand is placed on the bus. A bus enforces a fixed ordering on allcommands placed on it. This order is agreed upon by all devices in thesystem since they all observe the same commands. The devices can alsoagree, without special effort, on the final effect of a sequence ofcommands. This is a major advantage for a single-bus-basedmultiprocessor.

A single-common-bus design, however, limits the size of the systemunless one opts for lower system performance. The limits of technologytypically allow only a few devices to be connected on the bus withoutcompromising the speed at which the bus switches and, therefore, thespeed at which the system runs. If more master devices, such asprocessors and I/O agents, are placed on the bus, the bus must switch atslower speeds, which lowers its available bandwidth. Lower bandwidth mayincrease queuing delays, which result in lowering the utilization ofprocessors and lowering the system performance.

Another serious shortcoming in a single-bus system is the availabilityof a single data path for transfer of data. This further aggravatesqueuing delays and contributes to lowering of system performance.Although a single-system-bus design is the current design choice ofpreference for implementing coherence protocol, it cannot be employedfor a large-way SMP with many processors.

Once a decision is made to design a large-way, distributedmultiprocessor system with multiple buses, there are several designchallenges for ensuring efficient data transfers. The number ofconnections to centralized control units can become substantial. Pincount on the physical components becomes a significant limitation,especially in a system that supports a large address space with largedata transfers. Hence, it is generally desirable to limit the number ofsignals so as to limit the number of physically separate pins. Inaddition, an effort should be made to increase the efficiency of busarbitration and data transfers so as to decrease the number of deadcycles on the bus.

Therefore, it would be advantageous to have a large-way SMP design usingbus-based cache-coherence protocols with efficient bus utilization anddata transfers.

SUMMARY OF THE INVENTION

A distributed system structure for a large-way, symmetric multiprocessorsystem using a bus-based cache-coherence protocol is provided. Thedistributed system structure contains an address switch, multiple memorysubsystems, and multiple master devices, either processors, I/O agents,or coherent memory adapters, organized into a set of nodes supported bya node controller. The node controller receives transactions from amaster device, communicates with a master device as another masterdevice or as a slave device, and queues transactions received from amaster device. Since the achievement of coherency is distributed in timeand space, the node controller helps to maintain cache coherency. Inaddition, a bus arbiter in the node controller parks a data bus towardsa memory subsystem. The node controller does not use data bufferreservations. The data bus grant line to the memory controller isoverloaded to use it as a back-pressure, get-off-the-bus signal as wellas a normal data bus grant line. The fairness of the bus is therebyincreased by creating a mechanism for getting a “parked” device off thedata bus without the use of another dedicated signal between physicalcomponents. To ensure that the node controller may stream data to thememory subsystem, the bus is not reparked towards the memory subsystemuntil a configurable number of cycles after the data bus has beengranted to the node controller.

BRIEF DESCRIPTION OF THE DRAWINGS

The novel features believed characteristic of the invention are setforth in the appended claims. The invention itself, however, as well asa preferred mode of use, further objectives and advantages thereof, willbest be understood by reference to the following detailed description ofan illustrative embodiment when read in conjunction with theaccompanying drawings, wherein:

FIG. 1 is a block diagram depicting the basic structure of aconventional multiprocessor computer system;

FIG. 2 is a block diagram depicting a typical architecture;

FIG. 3 is a block diagram depicting an SMP computer system with threeprocessing units;

FIG. 4 is a block diagram depicting a distributed system structure for asymmetric multiprocessor system with supporting bus-basedcache-coherence protocol from the perspective of address paths withinthe SMP system;

FIG. 5 is a block diagram depicting a distributed system structure for asymmetric multiprocessor system with supporting bus-basedcache-coherence protocol from the perspective of data paths within theSMP system;

FIG. 6 is a block diagram depicting the address paths internal to a nodecontroller;

FIG. 7 is a diagram depicting the internal address paths of an addressswitch connecting node controllers and memory subsystems;

FIG. 8 is a diagram depicting a memory subsystem connected to theaddress switch of the distributed system of the present invention;

FIGS. 9A-9B are block diagrams depicting the data paths internal to anode controller;

FIGS. 10A-10B are block diagrams depicting the system structure fordetermining bus response signals for a distributed system structure;

FIGS. 10C-10D are block diagrams depicting the components whose signalsparticipate in the local and global cycles;

FIG. 11 is a state diagram showing the states that may be used in astate machine for parking a data bus towards the memory controller asviewed by a processor;

FIG. 12 is a flowchart describing a process of parking a data bustowards the memory controller as viewed from the perspective of thememory controller; and

FIG. 13 is a state diagram that is similar to the state diagram shown inFIG. 11 but modified to introduce a delay parking condition.

DETAILED DESCRIPTION OF THE PREFERRED EMBODIMENT

With reference now to FIG. 1, the basic structure of a conventionalmultiprocessor computer system 110 is depicted. Computer system 110 hasseveral processing units 112 a, 112 b, and 112 c which are connected tovarious peripheral devices, including input/output (I/O) agents 114,which accept data from and provide data to a monitor adapter 102 anddisplay monitor 105, keyboard adapter 104 and keyboard 107, and diskadapter 103 and permanent storage device 106, memory device 116 (such asdynamic random access memory or DRAM) that is used by the processingunits to carry out program instructions, and firmware 118 whose primarypurpose is to seek out and load an operating system from one of theperipherals (usually the permanent memory device) whenever the computeris first turned on. Processing units 112 a-112 c communicate with theperipheral devices by various means, including a bus 120. Computersystem 110 may have many additional components which are not shown, suchas serial and parallel ports for connection to peripheral devices, suchas modems or printers. Those skilled in the art will further appreciatethat there are other components that might be used in conjunction withthose shown in the block diagram of FIG. 1; for example, a displayadapter might be used to control a video display monitor, a memorycontroller can be used to access memory 116, etc. In addition, computersystem 110 may be configured with more or fewer processors.

In a symmetric multiprocessor (SMP) computer, all of the processingunits 112 a-112 c are generally identical; that is, they all use acommon set or subset of instructions and protocols to operate andgenerally have the same architecture.

With reference now to FIG. 2, a typical organization is depicted. Aprocessing unit 112 includes a processor 122 having a plurality ofregisters and execution units, which carry out program instructions inorder to operate the computer. The processor can also have caches, suchas an instruction cache 124 and a data cache 126. These caches arereferred to as “on-board” when they are integrally packaged with theprocessor's registers and execution units. Caches are commonly used totemporarily store values that might be repeatedly accessed by aprocessor, in order to speed up processing by avoiding the longer stepof loading the values from memory, such as memory 116 shown in FIG. 1.

Processing unit 112 can include additional caches, such as cache 128.Cache 128 is referred to as a level 2 (L2 ) cache since it supports theon-board (level 1) caches 124 and 126. In other words, cache 128 acts asan intermediary between memory 116 and the on-board caches, and canstore a much larger amount of information (instructions and data) thanthe on-board caches, although at a longer access penalty. For example,cache 128 may be a chip having a storage capacity of 256 or 512kilobytes, while the processor 112 may be an IBM PowerPC™ 604-seriesprocessor having on-board caches with 64 kilobytes of total storage.Cache 128 is connected to bus 120, and all loading of information frommemory 116 into processor 112 must come through cache 128. Although FIG.2 depicts only a two-level cache hierarchy, multi-level cachehierarchies can be provided where there are many levels of seriallyconnected caches.

In an SMP computer, it is important to provide a coherent memory system,that is, to cause writes to each individual memory location to beserialized in some order for all processors. For example, assume alocation in memory is modified by a sequence of writes to take on thevalues 1, 2, 3, 4. In a cache-coherent system, all processors willobserve the writes to a given location to take place in the order shown.However, it is possible for a processing element to miss a write to thememory location. A given processing element reading the memory locationcould see the sequence 1, 3, 4, missing the update to the value 2. Asystem that ensures that each processor obtains valid data order is saidto be “coherent.” It is important to note that virtually all coherencyprotocols operate only to the granularity of the size of a cache block.That is to say, the coherency protocol controls the movement of thewrite permissions for data on a cache block basis and not separately foreach individual memory location.

There are a number of protocols and techniques for achieving cachecoherence that are known to those skilled in the art. At the heart ofall these mechanisms for maintaining coherency is the requirement thatthe protocols allow only one processor to have a “permission” thatallows a write to a given memory location (cache block) at any givenpoint in time. As a consequence of this requirement, whenever aprocessing element attempts to write to a memory location, it must firstinform all other processing elements of its desire to write the locationand receive permission from all other processing elements to perform thewrite command. The key issue is that all other processors in the systemmust be informed of the write command by the initiating processor beforethe write occurs. To further illustrate how cache coherence isimplemented in multi-level hierarchies, consider FIG. 3.

With reference now to FIG. 3, an SMP computer system is depicted withthree processing units (140, 141, 142) consisting of processors (140 a,141 a, 142 a) each having an L1 cache (140 b, 141 b, 142 b), and L2cache (140 c, 141 c, 142 c), and finally, an L3 cache (140 d, 141 d, 142d). In this hierarchy, each lower-level cache (i.e., an L3 cache is“lower” than an L2 ) is typically larger in size and has a longer accesstime than the next higher-level cache. Furthermore, it is common,although not absolutely required, that the lower-level caches containcopies of all blocks present in the higher-level caches. For example, ifa block is present in the L2 cache of a given processing unit, thatimplies the L3 cache for that processing unit also has a (potentiallystale) copy of the block. Furthermore, if a block is present in the L1cache of a given processing unit, it is also present in the L2 and L3caches of that processing unit. This property is known as inclusion andis well-known to those skilled in the art. Henceforth, it is assumedthat the principle of inclusion applies to the cache related to thepresent invention.

To implement cache coherency in a system such as is shown in FIG. 3, theprocessors communicate over a common generalized interconnect (143). Theprocessors pass messages over the interconnect indicating their desireto read or write memory locations. When an operation is placed on theinterconnect, all of the other processors “snoop” this operation anddecide if the state of their caches can allow the requested operation toproceed and, if so, under what conditions. This communication isnecessary because, in systems with caches, the most recent valid copy ofa given block of memory may have moved from the system memory 144 to oneor more of the caches in the system. If a processor (say 140 a) attemptsto access a memory location not present within its cache hierarchy (140b, 140 c and 140 d), the correct version of the block, which containsthe actual value for the memory location, may either be in the systemmemory 144 or in one of the caches in processing units 141 and 142. Ifthe correct version is in one of the other caches in the system, it isnecessary to obtain the correct value from the cache in the systeminstead of system memory.

For example, consider a processor, say 140 a, attempting to read alocation in memory. It first polls its own L1 cache (140 b). If theblock is not present in the L1 cache (140 b), the request is forwardedto the L2 cache (140 c). If the block is not present in the L2 cache,the request is forwarded on to the L3 cache (140 d). If the block is notpresent in the L3 cache (140 d), the request is then presented on thegeneralized interconnect (143) to be serviced. Once an operation hasbeen placed on the generalized interconnect, all other processing units“snoop” the operation and determine if the block is present in theircaches. If a given processing unit, say 142, has the block of datarequested by processing unit 140 in its L1 cache (142 a), and the datais modified, by the principle of inclusion, the L2 cache (142 c) and theL3 cache (142 d) also have copies of the block. Therefore, when the L3cache (142 d) of processing unit 142 snoops the read operation, it willdetermine that the block requested is present and modified in the L3cache (142 d). When this occurs, the L3 cache (142 d) may place amessage on the generalized interconnect informing processing unit 140that it must “retry” its operation again at a later time because themost recently updated value of the memory location for the readoperation is in the L3 cache (142 d), which is outside of main memory144, and actions must be taken to make it available to service the readrequest of processing unit 140.

The L3 cache (142 d) may begin a process to push the modified data fromthe L3 cache to main memory 144. The most recently updated value for thememory location has then been made available to the other processors.

Alternatively, in a process called “intervention,” the L3 cache (142 d)may send the most recently updated value for the memory locationdirectly to processing unit 140, which requested it. The L3 cache maythen begin a process to push the modified data from the L3 cache to mainmemory. Processing unit 140, specifically its L3 cache (140 d),eventually represents the read request on the generalized interconnect.At this point, however, the modified data has been retrieved from the L1cache of processing unit 142 and the read request from processor 140will be satisfied. The scenario just described is commonly referred toas a “snoop push.” A read request is snooped on the generalizedinterconnect which causes processing unit 142 to “push” the block to thebottom of the hierarchy to satisfy the read request made by processingunit 140.

The key point to note is that, when a processor wishes to read or writea block, it must communicate that desire with the other processing unitsin the system in order to maintain cache coherence. To achieve this, thecache-coherence protocol associates, with each block in each level ofthe cache hierarchy, a status indicator indicating the current “state”of the block. The state information is used to allow certainoptimizations in the coherency protocol that reduce message traffic ongeneralized interconnect 143 and inter-cache connections 140 x, 140 y,141 x, 141 y, 142 x, 142 y. As one example of this mechanism, when aprocessing unit executes a read, it receives a message indicatingwhether or not the read must be retried later. If the read operation isnot retried, the message usually also includes information allowing theprocessing unit to determine if any other processing unit also has astill active copy of the block (this is accomplished by having the otherlowest-level caches give a “shared” or “not shared” indication for anyread they do not retry).

In this manner, a processing unit can determine whether any otherprocessor in the system has a copy of the block. If no other processingunit has an active copy of the block, the reading processing unit marksthe state of the block as “exclusive.” If a block is marked exclusive,it is permissible to allow the processing unit to later write the blockwithout first communicating with other processing units in the systembecause no other processing unit has a copy of the block. Therefore, ingeneral, it is possible for a processor to read or write a locationwithout first communicating this intention onto the interconnection.However, this only occurs in cases where the coherency protocol hasensured that no other processor has an interest in the block. Severaldetails of the exact workings of a multi-level cache coherence protocolhave been omitted in this discussion to simplify it. However, theessential aspects that bear on the invention have been described. Thoseaspects that bear on the invention have been described. Those aspectsnot described are well-known to those skilled in the art.

Another aspect of multi-level cache structures relevant to the inventionare the operations known as deallocations. The blocks in any cache aredivided into groups of blocks called “sets”. A set is the collection ofblocks in which a given memory block can reside. For any given memoryblock, there is a unique set in the cache that the block can be mappedinto, according to preset mapping functions. The number of blocks in aset is referred to as the associativity of the cache (e.g., 2-way setassociative means that, for any given memory block, there are two blocksin the cache that the memory block can be mapped into). However, severaldifferent blocks in main memory can be mapped to any given set.

When all of the blocks in a set for a given cache are full and thatcache receives a request, whether a read or write, to a memory locationthat maps into the full set, the cache must “deallocate” one of theblocks currently in the set. The cache chooses a block to be evicted byone of a number of means known to those skilled in the art (leastrecently used (LRU), random, pseudo-LRU, etc.). If the data in thechosen block is modified, that data is written to the next lowest levelin the memory hierarchy, which may be another cache (in the case of theL1 or L2 cache) or main memory (in the case of an L3 cache). Note that,by the principle of inclusion, the lower level of the hierarchy willalready have a block available to hold the written modified data.However, if the data in the chosen block is not modified, the block issimply abandoned and not written to the next lowest level in thehierarchy. This process of removing a block from one level of thehierarchy is known as an “eviction.” At the end of this process, thecache no longer holds a copy of the evicted block and no longer activelyparticipates in the coherency protocol for the evicted block because,when the cache snoops an operation (either on generalized interconnect143 or inter-cache connections 140 x, 141 x, 142 x, 140 y, 141 y, 142y), the block will not be found in the cache.

The present invention discloses a distributed hardware structure toovercome the limitations of a single common bus in a multiprocessorsystem while utilizing the properties of the single bus so that it doesnot require a modification to the bus protocol. The resulting system hasa scalable system size without compromising the mechanism of a knownsystem bus. The present invention is able to connect together a largenumber of devices in an SMP system and overcome the limitations of asingle-bus-based design.

Although the following description describes the invention with respectto the 6XX bus architecture, the present invention is not intended to belimited to a particular bus architecture as the system presented belowcan be applied to other bus architectures.

System Address Path Topology

With reference now to FIG. 4, a block diagram depicts a distributedsystem structure for a symmetric multiprocessor system with supportingbus-based cache-coherence protocol from the perspective of address pathswithin the SMP system. FIG. 4 displays a number of master devices thatcan initiate a command, such as a memory transaction. These masterdevices, such as processors, I/O agents, and coherent memory adapters,are distributed in clusters among a number of N groups called nodes.Each node is headed by a node controller into which its masters connect.

FIG. 4 shows nodes 410 and 420, which contain groupings of systemelements. The number of nodes may vary based on the configuration of thesystem. Node 410, also labeled as Node₀, contains processors 411 and412, also labeled as Processor P₀ and Processor P_(P−1), which are themasters for Node 410. Each node controller has multiple standardbidirectional processor address-data buses over which masters areconnected into the distributed system. Processors 411 and 412 connect tonode controller 415, also labeled as Node Controller NC₀, via buses 413and 414, also labeled as P₀Bus and P_(P−1)Bus, respectively. Node 420,also labeled as Node_(N−1), contains processor 421 and I/O agent 422,which are the masters for Node 420. Processor 421 and I/O device 422connect to node controller 425, also labeled as Node Controller NC_(N−1)via buses 423 and 424, respectively. The number of masters per node mayvary depending upon the configuration of the system, and the number ofmasters at each node is not required to be uniform across all of thenodes in the system.

The node controller constitutes the physical interface between a masterand the rest of the system, and each node controller in the systemcontains all of the necessary logic to arbitrate for individualprocessor buses and to communicate with its local masters as anothermaster or as a slave, i.e. a device that accepts master commands andexecutes them but does not generate master commands. A processor sends acommand into the system via its local node controller. Although FIG. 4shows one master per port, multiple masters per port are possible givenan appropriate arbitration scheme on the bus of that port. For example,processor 411 could be one of many processors connected to bus 413.However, if more processors are connected to a single port, then theiraddress bus will perform more slowly in terms of bus cycle time.

Alternatively, one of the masters of Node 420 may include a coherentmemory adapter that provides communication with another data processingsystem that maintains cache coherence. The coherent memory adapter maybe proximate or remote and may occupy a port of a node controller tosend and receive memory transactions in order to behave as amaster/slave device in a manner similar to an I/O agent. As one example,another node controller from another data processing system may also beconnected to the coherent memory adapter so that data processing systemsthat employ the present invention may be chained together.

Node controllers 415 and 425 are connected to a device called an addressswitch (ASX) via pairs of unidirectional address-only buses. Buses 416and 417, also labeled AOut₀ and AIn₀, respectively, connect nodecontroller 415 to address switch 430. Buses 426 and 427, also labeledAOut_(N−1) and AIn_(N−1), respectively, connect node controller 425 toaddress switch 430. As shown, buses AOut_(X) carry addresses from thenode controllers to the address switch, and buses AIn_(X) carryaddresses from the address switch to the node controllers.

Address switch 430 has additional unidirectional address bus connections431 and 432, also labeled as AIn_(N) and AIn_((N+S−1)), to memorycontrollers or memory subsystems 442 and 444, also labeled as memorysubsystem MS₀ and MS_(S−1). The memory controllers are assumed to beslave devices and have no ability to issue commands into the distributedsystem. The number of memory subsystems may vary depending upon theconfiguration of the system.

System Data Path Topology

With reference now to FIG. 5, a block diagram depicts a distributedsystem structure for a symmetric multiprocessor system with supportingbus-based cache-coherence protocol from the perspective of data pathswithin the SMP system. In a manner similar to FIG. 4, FIG. 5 displays anumber of master devices. These master devices are distributed inclusters among a number of N groups called nodes. Each node is headed bya node controller into which its masters connect. FIG. 5 shows nodes 510and 520 containing processors 511 and 512. Processors 511 and 512connect to node controller 515 via buses 513 and 514. Node 520, alsolabeled as Node_(N−1), contains processor 521 and I/O device 522 thatconnect to node controller 525, also labeled as Node Controller NC_(N−1)via buses 523 and 524, respectively.

The node controllers shown in FIG. 4 and FIG. 5 could be physically thesame system component but are described from different perspectives toshow different functionality performed by the node controllers. WhereasFIG. 4 shows address paths within the SMP system, FIG. 5 shows the datapaths within the SMP system. Alternatively, in a preferred embodiment,the address paths and data paths may be implemented with supportingfunctionality in physically separate components, chips, or circuitry,such as a node data controller or a node address controller. The choiceof implementing a node controller with separate or combined data andaddress functionality may depend upon parameters of other systemcomponents. For example, if the sizes of the buses supported within thesystem are small enough, both address and data functionality may beplaced within a single node controller component. However, if the busessupport 128 bits of data, then pin limitations may physically requirethe address and data functionality to be placed within separate nodecontroller components.

Alternatively, a separate node data controller may be further separatedinto multiple node data controllers per node so that each node datacontroller provides support for a portion of the node's data path. Inthis manner, the node's data path is sliced across more than one nodedata controller.

In FIG. 5, each node controller is shown connected to a plurality ofmemory controllers, such as memory subsystems MS₀ and MS_(S−1). Althougheach node controller is shown to connect to each memory controller viaan independent data bus, multiple nodes and/or multiple memorycontrollers may be connected on the same data bus if an appropriatearbitration mechanism is included. As with connecting a plurality ofmaster devices to a single node controller via a single bus, theswitching rate will be a function of the number of devices connected tothe bus. Node controller 515 connects to memory subsystem 542 via databus 516, and to memory subsystem 544 via bus 517, also labeled as N₀D₀and N₀D_(S−1), respectively. Node controller 525 connects to memorysubsystem 544 via data bus 527, and to memory subsystem 542 via data bus526, also labeled as N_(N−1)D_(S−1) and N_(N−1)D₀, respectively.

Instead of a single data bus that transfers data belonging to all of themasters, there are multiple data buses, each of which carries only asmall portion of the data traffic that would be carried if the masterswere connected to a single bus. In so doing, the component interfacesmay be clocked faster than would be possible with a single bus. Thisconfiguration permits the allocation of more data bus bandwidth permaster than would be possible on a single bus, leading to lower queueingdelays.

Node Controller Internal Address Paths

With reference now to FIG. 6, a block diagram depicts the address pathsinternal to a node controller. Node controller 600, also labeled NC_(X),is similar to node controllers 415 and 425 in FIG. 4 or node controllers515 and 525 in FIG. 5. Individual ports of node controller 600 havetheir own queues to buffer commands from masters as the commands enterthe node controller. A command may incur non-deterministic delay whilewaiting in these buffers for progressive selection toward the addressswitch.

Node controller 600 has bidirectional buses 601-604 that connect tomaster devices. Buses 601-604 connect to input boundary latches 609-612and output boundary latches 613-616 via bus transceivers 605-608. Inputboundary latches 609-612 feed buffers 617-620 that hold the commandsfrom the master devices. A command from a master device may consist of atransaction tag, transaction type, target or source address, and otherpossible related information. Buffers 617-620 may hold all informationrelated to a command, if necessary, or may alternatively hold only theinformation necessary for the functioning of the address path within thenode controller. The information held by the input buffers may varydepending on alternative configurations of a node controller. Buffers617-620 feed control unit/multiplexer 621 that selects one command at atime to send to the address switch via latch 622, transmitter 623, andbus 624, also labeled AOut_(X).

Node controller 600 receives commands from masters via buses 601-604 foreventual transmittal through boundary latch 622 and transmitter 623 tothe address switch via bus 624, also labeled bus AOut_(X). In acorresponding manner, node controller 600 accepts commands from theaddress switch via bus 625, also labeled bus AIn_(X), and receiver 626for capture in boundary latch 627, also labeled as FROM_ASX_BL. Thesecommands follow an address path through a fixed number of latches thathave a fixed delay, such as intermediate latch 628 and output boundarylatches 613-616, before reaching buses 601-604. In addition, commands tomaster devices also pass through a multiplexer per port, such as controlunits/multiplexers 629-632, that also have a fixed delay. In thismanner, commands arriving via bus 625 traverse a path with a fixed delayof a deterministic number of cycles along the path. In other words, afixed period of time occurs between the point when a command reacheslatch FROM_ASX_BL to the point at which each master device, such as aset of processors connected to the node controller, is presented withthe arriving command.

The arbiters for the ports connected to the masters are designed to givehighest priority to the node controllers driving the port buses. If amaster makes a request to drive a bus at the same time that the nodecontroller expects to drive it, the node controller is given highestpriority. In a preferred embodiment, to assist with this arbitrationscenario, a signal called “SnoopValid” (not shown) is asserted by theaddress switch ahead of the command being sent by the address switch.This allows the arbitration for the bus accesses between a nodecontroller and its masters to be completed early enough to ensure that acommand arriving from the address switch via the AIn_(X) bus does notstall for even one cycle while inside the node controller. Thisguarantees that the time period for the fixed number of latches alongthe AIn_(X)-to-P_(X)Bus paths actually resolve to a deterministic numberof cycles.

Control logic unit 633 is also presented with the incoming commandlatched into the FROM_ASX_BL latch for appropriate determination ofcontrol signals to other units or components within node controller 600.For example, control logic unit 633 communicates with buffers 617-620via control signals 634, control unit/multiplexer 621 via controlsignals 636, and control units/multiplexers 629-632 via control signals635 to select commands, resolve collisions, and modify fields ofcommands, including a command's type if necessary, in order to ensurethe continuous flow of commands within node controller 600. Controllogic unit 633 also receives other control signals 637, as appropriate.

Address Switch Internal Address Paths

With reference now to FIG. 7, a diagram depicts the internal addresspaths of an address switch connecting node controllers and memorysubsystems. Address switch 700 connects a set of four node controllersand two memory subsystems. Commands arrive at first-in first-out (FIFO)queues 721-724 from buses 701-704, also labeled AOut₀-AOut₃, viareceivers 709-712 and input boundary latches 713-716. These commands mayreside within a FIFO before being selected by control unit/multiplexer725. A command may experience a finite but non-deterministic number ofcycles of delays while sitting in the FIFO. Control logic unit 726 maycommunicate with control unit/multiplexer 725 and FIFOs 721-724 in orderto determine the selection of incoming commands. Control logic unit 726also receives other control signals 733, as appropriate.

Control unit/multiplexer 725 selects one command at a time to bebroadcast to the node controllers and memory subsystems over paths thatare deterministic in terms of the number of cycles of delay. In theexample shown in FIG. 7, commands are sent to the memory subsystems viaunidirectional buses 731 and 732, also labeled as buses AIn₄ and AIn₅,through output boundary latches 727 and 728 and transmitters 729 and730. Commands are sent to node controllers via unidirectional buses705-708, also labeled as buses AIn₀-AIn₃, through output boundarylatches 717-720 and transmitters 741-744. In this example, there is onlya single cycle of delay at the output boundary latches 717-720, 727, and728.

From the descriptions above for FIGS. 4-7, it may be understood that atransaction is issued by a master device via its bus and port to itsnode controller. The node controller will provide some type of immediateresponse to the master device via the bus and may queue the transactionfor subsequent issuance to the rest of the system. Once the transactionis issued to the rest of the system, the address switch ensures that thetransaction can be broadcast to the rest of the system with a knownpropagation delay so that the other devices may snoop the transaction.

According to the distributed system structure of the present invention,each of the devices within the system would be able to see thetransaction in the same cycle and provide a coherence response withinthe same cycle. The address switch is able to broadcast a transaction toall node controllers, including the node controller of the nodecontaining the device that issued the transaction. Appropriate logic isembedded within each node controller so that a node controller maydetermine whether the incoming transaction being snooped was originallyissued by a device on one of its ports. If so, then the node controllerensures that the bus on the port that issued the transaction is notsnooped with a transaction that was received from that port. Otherwise,the device may get “confused” by being snooped with its own transaction.If the device were to receive a snoop of its own transaction, then thedevice may issue a response indicating a collision with its originaltransaction. If that were the case, since the original transaction isactually the transaction that is being snooped, then the “collision”would never be resolved, and the transaction would never complete.

More details of the manner in which the transactions are issued andcompleted are provided below.

Memory Subsystem Internal Address Paths

With reference now to FIG. 8, a diagram depicts a memory subsystemconnected to the address switch of the distributed system of the presentinvention. FIG. 8 shows memory subsystem 800, also labeled memorysubsystem MS_(X). Memory controller 801 within memory subsystem 800receives a command from the address switch via unidirectional bus 802,also labeled as bus AIn_(X), through a number of latches FD 803, whichis merely a fixed delay pipe. In this manner, a command sent by theaddress switch experiences a fixed number of cycles of delay before thecommand is made available to the memory controller.

As shown previously, a command arriving at a node controller via busAIn_(X) traverses a deterministic delay path from its capture in theFROM_ASX_BL latch to its presentation to a master device. In a similarmanner, a command traverses a deterministic delay path from the controlunit/multiplexer within the address switch to the fixed delay pipewithin the memory subsystem. If the delay of the latches FD 803 withinthe memory subsystem is adjusted to the appropriate value, it can beensured that the memory controller is presented with a command at thesame time that the masters connected to the ports of the nodecontrollers are presented with the same command. Hence, there is adeterministic number of cycles between the point at which the controlunit/multiplexer within the address switch broadcasts a transaction andthe point at which the masters and memory controllers receive thecommand.

Since only a small number of masters are connected to each port of anode controller, the speed at which each bus is connected to these portsmay be operated is independent of the total number of ports in thesystem. For example, if a single master is connected to each port, itsbus can be run in point-to-point mode at the best possible speed. Hence,the distributed structure of the present invention is able to scalewell-understood and easier-to-verify bus-based cache-coherent protocolsfor multiprocessors to enhance the bandwidth of the system.

Node Controller Internal Data Paths

With reference now to FIGS. 9A-9B, block diagrams depict the data pathsinternal to a node controller. Node controller 900, also labeled NC_(X),is similar to node controllers 415 and 425 in FIG. 4 or node controllers515 and 525 in FIG. 5. Individual ports of node controller 900 havetheir own queues to buffer data from masters as data enters the nodecontroller. Data may incur non-deterministic delay while waiting inthese buffers for progressive movement toward destinations.

Node controller 900 has bidirectional buses 901-904, also labeledP_(X)Bus, that connect to master devices. Buses 901-904 connect to inputboundary latches 909-912 and output boundary latches 913-916 via bustransceivers 905-908. Input boundary latches 909-912 feed data buffers917-920 that hold the data from the master devices.

Incoming data from one of the node controller's ports may be directed toa memory subsystem or another cache. In the examples shown in FIGS.9A-9B, which continues the example shown in FIG. 6, incoming data fromone of the node controller's ports may be directed to one of threelocations: memory subsystem MSo, memory subsystem MS_(S−1), or acache-to-cache FIFO (FIFO C2C) for forwarding data within the node. Withthe FIFO C2C mechanism, each node is able to transfer data from one ofits ports to another port, thereby allowing the transfer of data fromone master to another. Buffers 917-920 feed multiplexers 925-927 thatselect a data source for forwarding data. Control logic unit 939provides control signals for multiplexer 925 to select data to be sentto memory subsystem MSo and for multiplexer 926 to select data to besent to memory subsystem MS_(S−1). Node controller 900 sends data frommultiplexers 925 and 926 through boundary latches 931 and 933 andtransceivers 935 and 936 to memory subsystem MS₀ and memory subsystemMS_(S−1), via bidirectional buses 937 and 938, also labeled N_(X)D₀ andN_(X)D_(S−1). Control logic unit 939 provides control signals formultiplexer 927 to select data to be forwarded within the node. Data isthen queued into FIFO 928.

In a corresponding manner, node controller 900 accepts data throughtransceivers 935 and 936 and boundary latches 932 and 934 from memorysubsystem MS₀ and memory subsystem MS_(S−1) via bidirectional buses 937and 938. Data is then queued into appropriate FIFOs 929 and 930. Datafrom FIFOs 928-930 pass through a multiplexer per port, such as controlunits/multiplexers 921-924. Control logic unit 939 provides controlsignals for multiplexers 921-924 to select data to be sent to the masterdevices. Control logic unit 939 also receives other control signals 940,as appropriate. Hence, the node controller has arbitration logic fordata buses and is self-sufficient in terms of controlling the datatransfers with parallelism. In this manner, the distributed systemstructure of the present invention is able to improve system datathroughput.

Response Combination Block (RCB)

With reference now to FIGS. 10A-10B, block diagrams depict the systemstructure for determining bus response signals for a distributed systemstructure similar to that shown in FIG. 4 and FIG. 5. FIG. 10A and FIG.10B show the connectivities of devices in the distributed systemstructure of the present invention with a control logic block forcombining bus signals (responses) AStat and AResp, respectively. For thesake of clarity, the AStat signals and the AResp signals have been shownseparately. It should again be noted that I/O agents may act as masterdevices connected to the ports of the node controllers shown in FIG. 10Aand FIG. 10B.

As shown in FIG. 10A, processors 1001-1004, also labeled P_(X), haveunidirectional AStatOut signals 1005-1008, also labeledP_(X)N_(X)AStOut, and AStatIn signals 1009-1012, also labeledP_(X)N_(X)AStIn, connecting the processors to Response Combination Block(RCB) 1000. The slave devices, such as memory subsystems 1005 and 1006,also labeled MS_(X), connect to the RCB with AStatOut signals 1013 and1014, also labeled M_(X) _(—) AStOut, and with AStatIn signals 1015 and1016, also labeled M_(X) _(—) AStIn. Node controllers 1017 and 1018,also labeled NC_(X), also connect to the RCB via a similar set of perport unidirectional AStatOut signals 1019-1022, also labeledN_(X)P_(X)AStOut, and AStatIn signals 1023-1026, also labeledN_(X)P_(X)AStIn. Address switch 1027, also labeled ASX, participates indetermining the proper logic for system processing of a transaction bysupplying broadcast signal 1028 and transaction source ID 1029, which isan encoding of a node identifier together with a port identifier withinthe node through which a master device issued a transaction to thesystem.

As shown in FIG. 10B, processors 1001-1004 have unidirectional ARespOutsignals 1055-1058, also labeled P_(X)N_(X)AReOut, and ARespIn signals1059-1062, also labeled P_(X)N_(X)AReIn, connecting the processors toRCB 1000. Memory subsystems 1005 and 1006 connect to the RCB withARespIn signals 1065 and 1066, also labeled M_(X) _(—) AReIn. Memorysubsystems 1005 and 1006 do not connect with ARespOut lines, which arenot driven by these slave devices. Node controllers 1017 and 1018 alsoconnect to the RCB via a similar set of per port unidirectional ARespOutsignals 1069-1072, also labeled N_(X)P_(X)AReOut, and ARespIn signals1073-1076, also labeled N_(X)P_(X)AReIn. Again, address switch 1027participates in determining the proper logic of a transaction bysupplying broadcast signal 1078 and transaction port ID 1079. RCB 1000supplies a HOLDTAG signal, such as signals 1091 and 1092, to each nodecontroller in certain circumstances, as explained further below.

As is apparent from FIGS. 10-10B, a set of AStatIn/AStatOut signals andARespIn/ARespOut signals to/from a master device is paired with asimilar set of AStatIn/AStatOut signals and ARespIn/ARespOut signalsto/from its node controller. This pairing is done on a per port basis.As discussed above, each port in the example is shown with a singlemaster device connected to each port. However, if more than one masterdevice were connected per port, then the pairs of AStatIn/AStatOutsignals and ARespIn/ARespOut signals are used by the set of masterdevices connected to the bus on that port as in a standard single busconfiguration.

In the preferred embodiment, RCB combines the AStatOuts and ARespOutsfrom various source devices and produces AStatIn and ARespIn signals perthe 6XX bus specification, as described in IBM Server Group Power PC MPSystem Bus Description, Version 5.3, herein incorporated by reference.The RCB receives the AStatOuts and ARespouts signals and returnsAStatIns and ARespIns, respectively. Not all of the devices receive thesame responses for a particular transaction. The signals received byeach device are determined on a per cycle basis as described in moredetail further below.

Local/Global cycles

During any given system cycle, a master device at a port may be issuinga transaction over its port's bus for receipt by its node controller orthe node controller may be presenting the master device with atransaction forwarded by the address switch in order to snoop thetransaction. When the master device is issuing a transaction, the cycleis labeled “local,” and when the node controller is presenting atransaction, the cycle is labeled “global.”

As described above, the address switch broadcasts one transaction at atime to all of the node controllers, and there is a fixed delay betweenthe time the address switch issues such a transaction and the time itappears at the ports of each node controller. Under this regime, after anode controller has received a broadcast transaction from the addressswitch and then, a predetermined number of cycles later, is presentingthe transaction to the devices on the buses of the ports of the nodecontroller during a cycle, all node controllers are performing the sameaction on all of their ports during the same cycle, except for oneexception, as explained below. Thus, when there is a global cycle beingexecuted on the bus of one of the ports, global cycles are beingexecuted on all the ports in the system. All remaining cycles are localcycles.

During local cycles, activity at a port is not correlated with activityat other ports within the system. Depending on whether or not a deviceneeded to issue a transaction, the local cycle would be occupied orwould be idle. Hence, a global cycle occurs when a transaction is beingsnooped by all the devices in the system, and only a local cycle may beused by a device to issue a transaction.

Operation of RCB During Local Vs Global Cycles

Given that the entire system's cycles are “colored” as either local orglobal, the response generation, the response combination, and theresponse reception cycles, which occur after a fixed number of cyclessubsequent to the issuance of a transaction, are similarly labeled localresponse windows or global response windows. For this reason, the RCB'sresponse combination function is correspondingly considered to be ineither local or global mode during a given cycle. During local cycles,the RCB combines responses on a per port basis. That is, the RCBcombines the response of a port and the response that the nodecontroller produces corresponding to that port. During global cycles,the RCB combines responses from all the ports and node controllers inthe system (again, except for one port, as explained below).

To achieve proper switching between local and global combination modes,the RCB is provided with a signal indicating the broadcast of atransaction by the address switch to the node controllers, shown asbroadcast signal 1028 in FIG. 10A, as well as the transaction source IDsignal 1029. Configuration information stored in the RCB indicates theexact cycle in which the combination of responses is to be performed forthe broadcast transaction after the arrival of the broadcast transactionsignal. In this manner, for each global cycle, the RCB is orchestratedto combine responses from appropriate sources.

Primary Vs Secondary Local cycles

A processor may issue a transaction only during local cycles. Forcertain types of transactions, the processor issues the transaction onlyonce. For certain other types of transactions, the processor might berequired to issue the transaction multiple times. The processor isdirected by its node controller, in conjunction with the RCB, throughthe use of the AStatIn/AStatOut signals and the ARespIn/ARespOut signalsas to the actions that should be performed.

The local cycles in which a processor issues transactions for the firsttime are labeled “primary local cycles” whereas all other local cyclesare labeled “secondary local cycles”. In the 6XX bus architecture, asecondary transaction is marked by the “R” bit being set to “1”. Inother words, its response-related cycles get labeled primary orsecondary in the proper manner corresponding to the transactionissuance.

Achievement of Coherence by Snooping in a Temporally and SpatiallyDistributed Manner

From the foregoing description, it should be obvious that processors anddevices see transactions from other processors and devices during cyclesdifferent than the cycle in which are issued to the system. This isunlike the situation with a snooping protocol in a single busenvironment in which all the devices in the system observe a transactionat the same time that-it is-issued and simultaneously produce acoherence response for it and in which the originator of the transactionreceives the response at that same time. Thus, in the current system,the achievement of coherence is both distributed in time and distributedin space, i.e. across multiple cycles and multiple buses connected tomultiple node controllers.

In using the distributed system structure, it is important to achieveglobal coherence in an efficient manner. To do so, all transactions aresorted into two categories: (1) transactions for which it is possible topredict the global coherence response and deliver it in the primaryresponse window; and (2) transactions for which it is necessary to snoopglobally before the ultimate coherence response can be computed.

In the first case, the node controller accepts the transaction andissues a global coherence response to the issuing entity in the primaryresponse window. The node controller then takes full responsibility ofcompleting the transaction in the system at a later time and achievingthe global response.

In the second case, the node controller takes three steps. First, thenode controller accepts the transaction and delivers a primary responsethat indicates postponement of achievement and delivery of the globalresponse. In the 6XX bus architecture, this response is the “Rerun”response. Second, at a subsequent time, the node controller achieves aglobal coherence response for that transaction. And third, the nodecontroller requests that the processor issue a secondary transaction anddelivers the global response in the secondary response window. In the6XX bus architecture, the request to the processor to issue a secondarytransaction is made by issuing it a Rerun command with a tagcorresponding to the original transaction. The processor may then usethe tag to identify which of its transactions should be rerun.

Rerun Commands and Secondary Responses

As noted above, a transaction accepted from a device is snooped to therest of the system. During such a snoop, the device that issued thetransaction is not snooped so that the device does not get confused bybeing snooped with its own transaction.

In fact, for transactions in the first case above, i.e. transactions inwhich the node controller accepts the transaction and issues a globalcoherence response to the issuing entity in the primary response window,the port corresponding to the device that issued the transaction is keptin the local mode in the transaction's snoop cycle so that the processormay issue another transaction. As stated above, during the responsewindow corresponding to the transaction's snoop cycle, the RCB isconfigured to combine responses from all sources other than the port onthe node controller that issued the transaction. The node controller isthen able to supply a primary or secondary response over that port ifthe processor chooses to issue a transaction.

For transactions in the second case above, i.e. transactions for whichit is necessary to snoop globally before the ultimate coherence responsecan be computed, the node controller keeps the particular port in localmode but issues it a Rerun transaction. The control unit/multiplexerfeeding the outgoing boundary latch at the port allows the nodecontroller to achieve this functionality.

Alternatively, the node controller may choose to not be as aggressive,and instead of letting the device issue a transaction, the nodecontroller might itself issue a null or rerun transaction, as required,to the device in the cycle during which the device's transaction isbeing snooped in the rest of the system.

With reference now to FIGS. 10C-10D, block diagrams depict thecomponents whose signals participate in the local and global cycles.FIG. 10C shows the signals which are considered by the RCB during aglobal cycle. In the example shown, the signals for a single masterdevice, processor 1001, do not participate in the determination by theRCB of the appropriate signals to the other devices, node controllers,and memory subsystems for the global response. The signals for processor1001 are paired with the corresponding signals from its node controller,which are also not considered for the global response. From theperspective of processor 1001, it is kept in a local cycle while atransaction issued by processor 1001 is snooped by the rest of thesystem. As noted earlier, although a processor is depicted, the signalsare considered on a per port basis, and the bus of a particular port iskept in a local cycle while the rest of the system is in a global cycle.

FIG. 10D shows the signals which are considered by the RCB during alocal cycle. In the example shown, the signals from a single masterdevice, processor 1001, participate in the determination by the RCB ofthe appropriate signals to be returned to processor 1001 and its nodecontroller. Signals from the other devices, node controllers, and memorysubsystems may be simultaneously participating in the response for theglobal response. The signals for processor 1001 are paired with thecorresponding signals from its node controller, which also do not affectthe global response. From the perspective of processor 1001, it mayissue another transaction while its other transaction is snooped by therest of the system. For the sake of clarity, signals from the addressswitch are not shown for the local cycle, although the RCB uses thesesignals to determine which port to place into the local cycle.

Achieving Correct Order Among Bus Memory Transactions

For a computer system to work correctly, certain memory accesstransactions and other types of transactions issued by master deviceshave to be ordered correctly and unambiguously. In a system with asingle system bus, this task is trivially achieved since the order inwhich the transactions are presented on the bus is the order imposed onthose transactions. However, in a distributed system with multiplebuses, the task demands that an order be imposed on the transactionsqueued throughout the system. The distributed architecture of thepresent invention allows a correct and unambiguous order to be imposedon a set of transactions. The invention also offers an efficient meansof achieving the order so that a snooping, hardware cache-coherenceprotocol can be supported.

When devices in an SMP system access memory, either under the influenceof programs or control sequences, they issue memory transactions. Thedevices may also issue other bus transactions to achieve coherence,ordering, interrupts, etc., in the system. These transactions canusually complete in parallel without interference from othertransactions. However, when two transactions refer to addresses withinthe same double word, for example, they are said to have “collided,”according to the 6XX bus terminology, and the two transactions must becompleted in some specific order. In some cases, either completion orderis acceptable, and at other times, the order is fixed and is implied bythe types of transactions. For instance, if a read transaction and aWrite transaction attempt to access an address declared as MemoryCoherence Not Required, any order of completion for the two transactionsis acceptable. However, if they refer to a cachable address to bemaintained coherent, the order of completion must appear to be the writefollowed by the read.

Means of Imposing a Default Order on Transactions

In the distributed SMP system described in FIGS. 4-10D, multipleprocessors and other devices can issue transactions simultaneously overthe multiple buses in the system. Thus, at the outset, there isambiguity regarding the order of the transactions as they are issued. Asthey flow through the system, as a first step, the system imposes a“heuristic order of arrival” over them that is reasonable and fair. Thispreliminary order is not necessarily the order in which the transactionseventually complete in the system. If two colliding transactions aresimultaneously active in the system, the one that ranked “earlier of thetwo” by the heuristic order of arrival will be slated to be completedfirst if coherence does not require otherwise.

As soon as commands enter the system, they are “registered” by the nodecontrollers, i.e. they are stored by the node controllers and areavailable for analysis and collision checks. Node controllers send oneof the registered transactions at a time to the address switch. Theaddress switch chooses one transaction at a time with a fair arbitrationamong the transactions sent to it and then broadcasts the chosentransaction back to the node controllers and to the memory subsystems.The address portion of the transaction broadcast by the address switchis first latched inside the node controller in the boundary latchFROM_ASX_BL. As described above, in any cycle, a unique transaction islatched in FROM_ASX_BL at all node controllers and memory subsystems,and all other registered transactions that have entered until that cycleand are still active, including the transaction currently inFROM_ASX_BL, can “see” this transaction. These two properties are usedto define the order of arrival of transactions using the followingreasonable and fair heuristic: the order of arrival of a transactioninto the system is the same as the order of its arrival at FROM_ASX_BL.

When a transaction arrives in FROM_ASX_BL for the first time, it ismarked as being “snooped,” to indicate the fact that in a fixed numberof cycles following the current cycle, the transaction will be presentedfor snooping, for the first time, to all the devices in the system. Thefollowing rule is used to assign a transaction its relative position inthe order of transactions to be completed, irrespective of the actualtime it entered the system: a registered transaction that already ismarked as snooped is nominally defined to have entered the systemearlier than the current transaction in FROM_ASX_BL. The ones that havenot been marked as snooped are nominally defined to have entered thesystem later than the current transaction in FROM_ASX_BL.

Method for Achieving the Correct Completion Sequence for Transactions

The transaction in FROM_ASX_BL stays there for one cycle. During thatcycle, the transaction is compared with every transaction currentlyregistered in the entire system for detection of collision and orderingdecision. There could be two sets of results of each of these pairwisecomparisons: one that affects the completion of the transactioncurrently in FROM_ASX_BL and the second that affects the completion ofsome other transaction.

Each comparison results in a decision to either allow the currentpresentation of the transaction in FROM_ASX_BL for snooping to complete,or to postpone its completion to a later time. The postponement iseffected via the computation of an AStat Retry signal or an AResp Retrysignal, as is appropriate. These signals from individual comparisons arecombined on a per node basis inside the node controller. A decision topostpone gets the highest priority, so even a single comparison callingfor postponement wins and results in the node voting to postpone thetransaction. Only if all comparisons-within a node vote to allow thecurrent snoop to complete does the node decide to let the transactioncomplete.

The combined AStat Retry and AResp Retry signals are encoded by the nodecontroller into the AStat Retry and ARespRetry codes and are submittedto the RCB for participation in the global AStat and AResp windows ofthe transaction being snooped. During these windows, responses from allthe devices, other than the device that issued the transaction, and nodecontrollers are combined by the RCB to produce a global response whichis returned to all the participants, as explained with respect to FIGS.10A-10D above. Again, at this global level, a retry response has thehighest priority (barring an error code) and will be the final responseif any of the input responses was a retry. The effect of a global retryresponse is cancellation of the current snoop of the transaction. Uponsensing a global retry response for the transaction, the node controllerin which the transaction is registered either reissues the transactionfor global snoop or retires the original transaction from which the saidtransaction was derived.

These global retries can be repeated until the correct order isachieved.

If, for any reason, a transaction receives a retry response, its snoopedmarking is reset, and it thus loses its present nominal position in thetransaction order in the system. When it returns for snoop, thetransaction gets a new position, according to the rule above. Themechanism does not necessarily prohibit the possibility of the reissuedtransaction being ordered behind another transaction that entered thesystem after it. If, on the other hand, the current transactioncompletes, it may cause other transactions to get retried.

Generalizing Bus Arbitration and Protocols

As described above, the node controller constitutes the physicalinterface between a master and the rest of the system, and each nodecontroller in the system contains all of the necessary logic toarbitrate for individual processor buses and to communicate with itslocal masters as another master or as a slave, i.e. a device thataccepts master commands and executes them but does not generate mastercommands. A processor sends a command into the system via its local nodecontroller, which then queues the commands and assumes responsibilityfor completing the commands in some form.

The following sections describe operational modes of a bus between aprocessor and a node controller or between a node controller and thememory subsystem. However, the discussion of the bus protocols can begeneralized by noting that the operations of a node controller, at leastfrom the perspective of arbitrating for the bus, may be replaced by aprocessor with a similar arbiter. This observation both simplifies thefollowing discussions and generalizes the present invention asoperational modes between two processors or between a processor and amemory subsystem.

It should be noted that one of ordinary skill in the art wouldunderstand that a bus arbiter may be included within the processor or,alternatively, may be a physically separate component with connectionsto the processor.

Fair Data Bus Parking Protocol Without Data Buffer Reservations on theReceiving End

Under most bus implementations, there is normally a presumption thatdata transfers from a memory controller back to a processor have anassociated data buffer reserved for that data. Thus, if N requests fordata are sent to the memory controller, there are N data buffersreserved in the requester for the data when it becomes available. Totake advantage of this, data bus arbitration schemes usually allow bushogging by the memory controller. This means that once the memorycontroller is given a data bus grant, the memory controller may continueto send data as long as it has data to send.

The present invention preserves some of this ability to stream dataduring bus hogging while not requiring data buffer reservations. Therecan be more data requests outstanding than there are data buffers to putthe data. The present invention achieves this by using the data busgrant line to the memory controller as a back-pressure signal or a “GetOff The Bus” signal as well as a normal data bus grant line. It alsoincreases the fairness of such a bus by creating such a mechanism forgetting a “parked” device off the data bus without the use of anotherdedicated signal between chips. This allows the processor to get ontothe bus and send data to memory more quickly, increasing fairness of useof the data bus.

The protocol involves asserting the data bus grant line continuously tothe memory controller after a data bus request is received, i.e. thedata bus grant is “parked” towards the memory controller. In thiscontext, a parked bus is a bus in which bus grants are given for the buseven when not requested. The memory controller is then free to send asmuch data back-to-back, i.e. “brick-walled”, as the memory controllerneeds to free up its data buffers. However, on the receiving end in theprocessor, which also contains the arbiter, an internal buffer-fullsignal is sent to the arbiter by other logic in the processor when acritical buffer-full condition is reached. This may occur if the datacannot be moved out of these data buffers as fast as it is being putinto the buffers. The number of buffers remaining before reaching abuffer-full condition varies with the bus design and system timings, buta typical count may be one or two buffers remaining. When this bufferfull signal is asserted, the arbiter will remove the data bus grant fromthe memory controller, acting as a “Get Off The Bus” signal. The memorycontroller will then complete any committed transactions that cannot beaborted, release the data bus, and reassert the data bus request line ifit still has data to transfer. It should be noted that if the arbiter isexternal to the processor, a buffer-full signal would be requiredbetween the processor and the arbiter.

The grant is also removed when the processor requests the data bus tosend data to the memory controller. After the grant has been given tothe processor, the memory controller is “reparked”—in other words, thememory controller is immediately given a continuous grant again, whetheror not it has requested the data bus. This decreases the latency beforethe memory controller can send data again by removing the delaysassociated with the normal request/grant protocol. There is a penalty tothe processor's latency to get the data bus if the processor decides torequest the data bus again. The latency is introduced by the timerequired to “unpark” the data bus, i.e., remove the continuous grant andthen wait to see if the memory controller uses any of the grants justgiven. This penalty is considered acceptable in the interest of speedingup the return of data to the processor from the memory controller.

The memory controller must not consider any of the grants given duringthe continuous grant phase to be pending grants that may be used afterthe current grant expires if it already owns the bus. If, on the otherhand, the memory controller does not have a grant or ownership of thebus when it receives a grant, it may then consider the grant to bepending.

A state machine may be used to keep track of which device has the olderoutstanding request for the data bus to further ensure fairness of useof the data bus. In prior art implementations, the memory controller isgiven complete priority to send data back to the processor. This canlead to deadlocks in a multiprocessor system where processors are notdirectly attached to the memory controller but are instead attached vianode controllers which act on their behalf. This state machinecontributes to the prevention of such deadlocks.

A two-state machine keeps track of which request is older. It is setwhenever the processor requests the data bus at a time when either thememory controller is not already requesting the data bus or is currentlyreceiving a continuous data bus grant regardless of the state of thememory controller request line. It is cleared when the processorreceives a data bus grant.

The continuous grant may be achieved by adding the two-state statemachine to an otherwise normal data bus arbiter, which may have its ownstate machines. The two-state state machine is set when a grant is givento the memory controller and the processor is not requesting the databus and the data buffers are not full. The two-state state machine iscleared whenever the processor is requesting the data bus or the databuffers are full. When this bit is set, it causes the memorycontroller's data bus grant line to be asserted. However, it does notinfluence any state machines in the otherwise normal arbiter. To do sowould confuse those other state machines into thinking there are far toomany grants outstanding in a system that normally allows one activegrant and one pending grant but not more. Thus, this invention is easilyadded to a normal, well debugged, data bus arbiter design with nochanges to the heart of the original arbiter. The fulfillment of thedata bus parking protocol is explained in more detail further below bydescribing how the two-state state machine may be added to an otherwisenormal data bus arbiter.

The memory controller must adhere to the following algorithm whenrequesting the data bus and using data bus grants for this protocol towork properly and thus avoid data buffer overflow:

1. If the data bus is not in use at the time that the memory controllerdetermines that the data bus is needed, then the data bus request lineis asserted—this is a start-from-idle condition.

2. If the data bus is in use by the memory controller and if the databus grant to the memory controller is already asserted, then“brick-wall” the next data transfer against the current one.

3. If the data bus is in use by the memory controller and the data busgrant to the memory controller is not asserted, then assert the data busrequest line after the data bus has been released by the memorycontroller.

These conditions would require some modification to a memory controlleroriginally designed to work with a parked data bus that has data bufferreservations on the receiving (processor) end. Specifically, conditions2 and 3 would not normally be implemented. This invention overloads themeaning of the data bus grant line to be able to avoid adding signallines.

With reference now to FIG. 11, a state diagram shows the states that maybe used in a state machine for parking a data bus towards the memorycontroller as viewed by a processor. As noted above, the parked data busgrant may be achieved by adding two-state state machine 1100 to anotherwise normal data bus arbiter, which may have its own statemachines. In state 1102, the data bus is not parked. State 1102 may beviewed as overlapping a state in a normal bus arbiter in which thenormal bus arbiter has not given a data bus grant to the memorycontroller. Before state machine 1100 may become “active”, i.e. beforethe conditions for parking the data bus may be considered, the normalbus arbiter must have given a data bus grant to the memory controller.If the memory controller has not first requested and received a data busgrant, it would be improper to park the data bus towards the memorycontroller as the memory controller may be offline or otherwiseinactive. Hence, state machine 1100 cannot move out of state 1102 untila flag is set, such as Park_Enabled, that informs the data bus arbiterthat it is acceptable to park the bus, or until the data bus grant hasbeen given to the memory controller via other state machines orconsiderations in the normal bus arbiter. The Park_Enabled flag isdescribed in more detail further below.

In addition, a defeat mechanism may be employed to disable the parkingalgorithm if desired. A flag could be set, such as Config_Park_Disable,to inform the data bus arbiter whether it is acceptable or unacceptableto attempt to park the bus. This easily defeats the parking mechanism ifit is unnecessary or undesired for some reason, such as debugging someother portion of a system.

The data bus arbiter loops in state 1102 waiting for the followingcondition to be true: (1) Parking is already enabled, i.e. thePark_Enabled flag is true, or the data bus grant has been given to thememory controller via other state machines or considerations in thenormal bus arbiter, i.e. DBG_(MC) is true; AND (2) the processor doesnot have a buffer-full condition, i.e. has not asserted a Buffer_Fullsignal to the arbiter or Buffer_Full is false; and there is no data busrequest from the processor, i.e. DBR_(PROC) is false; and the data busmay be parked, i.e. Config_Park_Disable is false. Under that condition,state machine 1100 moves from state 1102, data bus not parked, to state1104, data bus parked.

Once state machine 1100 reaches state 1104, the Park_Enabled flag is setto true to inform the data parking mechanism whether state machine 1100is “active”, i.e. whether the conditions for parking the data bus shouldbe considered. In other words, this ensures that state machine 1100 canbe moved from an initialization state and allows the data bus arbiter tobegin parking the data bus if the proper conditions for doing so exist.Once the Park_Enabled flag is set to true, there is no harm incontinuing to set the flag when state machine 1000 returns to state1104. As long as the condition remains that allowed the data bus to beparked, state machine 1100 loops in state 1104, i.e. the data busremains parked.

In state 1104, as long as the processor does not assert a data busrequest, i.e. DBR_(PROC) is false, and the processor does not have abuffer-full condition, i.e. has not asserted a Buffer_Full signal orBuffer_Full is false, then the data bus stays parked towards the memorycontroller as the data bus is not needed by the processor.

However, if either of these conditions changes, the data bus isunparked. In other words, the data bus arbiter loops in state 1102waiting for either of the following conditions to be true: (1) if theprocessor has a buffer-full condition, i.e. has asserted a Buffer_Fullsignal to the arbiter or Buffer_Full is true; OR (2) there is a data busrequest from the processor, i.e. DBR_(PROC) is true. Under either ofthese conditions, state machine 1100 moves from state 1104, data busparked, to state 1104, data bus not parked.

After state machine 1100 returns to state 1102, the arbiter returns tonormal ping-ponging of data bus arbitration between the processor andthe memory controller.

In this manner, state machine 1100 describes an algorithm for parking adata bus from the perspective of the processor in which the memorycontroller receives a data bus grant either during the normal course ofoperation of a data bus arbiter and its associated state machines or theparking of the data bus grant to the memory controller through the databus parking mechanism described above.

With reference now to FIG. 12, a flowchart describes a process ofparking a data bus towards the memory controller as viewed from theperspective of the memory controller. The process begins when the memorycontroller determines whether it has a new packet or next packet of datato transfer to the processor (step 1202). If not, then the memorycontroller loops back to step 1202, in which case it may besimultaneously performing other activities.

If there is a data packet to transfer, then the memory controllerdetermines whether the data bus busy signal is asserted (step 1204). Ifnot, then the memory controller asserts the data bus request signalimmediately (step 1206). The memory controller then determines whetherthe data bus grant is asserted (step 1208). If not, the memorycontroller loops back to step 1208 to wait for its assertion. If thedata bus grant is asserted, then the memory controller deasserts thedata bus request (step 1210) and begins transferring the first datapacket (step 1212). The process then loops back to step 1202 todetermine whether there is a next packet of data to be transferred. Inthis manner, steps 1206-1212 describe a standard operation of a datatransfer from a standard memory controller.

If the data bus busy signal is asserted in step 1204, then adetermination is made as to whether the data bus grant is asserted (step1214). If not, then the process loops back to step 1204 to wait for thedata bus grant. If so, then the memory controller can “brickwall” thenext data packet transfer without a pause between transfers (step 1216).Hence, step 1214 provides the indication from the arbiter to the memorycontroller with a “Get Off The Bus” signal in accordance with apreferred embodiment of the present invention.

In this manner, the memory controller looks at the data bus grant signalas a back-pressure signal from the processor to the memory controller.Thus, the data bus grant signal has meaningful content both on itstransition from “0” to “1” but also on its transition from “1” to “0”.This differs from a standard data bus grant signal that has significanceonly for its change from “0” to “1”. With the use of a standard data busgrant signal, memory controller would move from determining whether thedata bus busy signal was asserted in step 1204, and if so, would thenbrickwall the next data packet, i.e., the memory controller wouldnormally hog the data bus as the memory controller might expect theprocessor to have enough data buffer space reserved for the next datatransfer.

Increasing Performance of a Parked Data Bus in the Non-Parked Direction

As explained with respect to FIG. 11 and FIG. 12, a protocol wasdescribed to park a memory controller on a data bus connected to aprocessor that does not use data buffer reservations. In this protocol,the data bus is reparked toward the memory controller immediately afterthe processor receives a data bus grant. If the processor requests thedata bus again, it must incur a several cycle penalty before receivingits next grant to allow the arbiter to “age out” the parked grants tothe memory controller. Due to bus staging delays, it may take a fewcycles for the arbiter to see a data bus request from the memorycontroller, and it may take a few cycles to determine whether the memorycontroller has used the data bus request. Many data bus grants could be“stacked up” on the data bus to fulfill the single request from thememory controller, and the arbiter must wait a few cycles to see if thedata bus is actually used by the memory controller. In other words, thememory controller may ignore a pending data bus grant, and the arbiter,in essence, waits for the pending grant to time-out. This effectivelyprevents the processor from streaming data to the memory controller.

Instead, the reparking of the data bus towards the memory controller maybe delayed by a configurable number of cycles. This time delay allowsthe processor to request the bus again and receive a pending grantwithout delay if the next request from the processor closely follows thelast grant. This will allow the processor to send data with minimaladditional cycles between transfers.

By adding logic to a data bus arbiter between two entities, one being aprocessor or node controller, the other being a memory controller, thedata bus arbiter can be configured to give a continuous, or parked, databus grant to the memory controller after a configurable delay from thepoint at which the arbiter has given a grant to the processor, assumingthere are no other requests or grants outstanding in the system.

By adding a delay parking counter in the arbiter to count the number ofcycles since the processor last received a data bus grant, the arbiterwill not repark the data bus towards the memory controller as long asthe value in the delay parking counter is less than a configurablenumber of cycles. When the delay parking counter reaches theconfigurable delay value, the data bus will be reparked towards thememory controller. This delay allows another request to arrive from theprocessor and gives the processor the pending grant instead of thememory controller, thus removing a several cycle penalty that would beincurred by the processor so that the arbiter may determine if thememory controller will actually use the pending data bus grant once theprocessor releases the data bus. This allows the processor to send databack-to-back. The delay parking counter is reset to zero when theprocessor receives a grant to start the process all over again.

The delay parking counter has no effect if the memory controllerexplicitly requests the data bus. It will be reparked as soon aspossible without regard to the delay parking counter in that case.

With reference now to FIG. 13, a state diagram is shown that is similarto the state diagram shown in FIG. 11 but modified to introduce a delayparking condition. State machine 1300 in FIG. 13 is similar to statemachine 1100 in FIG. 11, and states 1302 and 1304 are similar to states1102 and 1104. However, state machine 1300 has an additional conditionbefore the data bus may be reparked towards the memory controller, asshown by the condition on Event B′. Each time that a data bus grant isgiven to the processor, i.e. DBG_(PROC) is true, an incrementer is resetto zero, and incremented after cycle thereafter. The arbiter isessentially blocked from examining the Parked_Enabled flag until thecounter reaches the configurable delay value. Thus, the processorreceives fairer treatment with respect to data bus grants.

Conclusion

The advantages of the present invention should be apparent in view ofthe detailed description provided above. The present invention allowsscaling of standardized and easier-to-verify bus-based cache-coherenceprotocols to a large-way, multiprocessor system whose large sizenormally would make physical buses inefficient media for communicationamong system components, such as processors, memory subsystems, and I/Oagents. By using the distributed system structure of the presentinvention, development of more complicated directory-based protocols,etc. are unnecessary. The present invention also allows componentinterfaces to be clocked faster than possible with a single bus, therebyenhancing the bandwidths of the component interfaces and resulting inhigher total system bandwidth and performance. The present inventionalso supports multiple data buses, thereby multiplying the databandwidth of the system and improving the efficiency of the processor.The data transfer parallelism of the present system also improves totalsystem data throughput.

As an additional advantage provided by the present system, a bus arbiterin the node controller parks a data bus towards a memory subsystem. Thenode controller does not use data buffer reservations. The data busgrant line to the memory controller is overloaded to use it as aback-pressure, get-off-the-bus signal as well as a normal data bus grantline. The fairness of the bus is thereby increased by creating amechanism for getting a “parked” device off the data bus without the useof another dedicated signal between physical components. To ensure thatthe node controller may stream data to the memory subsystem, the bus isnot reparked towards the memory subsystem until a configurable number ofcycles after the data bus has been granted to the node controller.

It is important to note that while the present invention has beendescribed in the context of a fully functioning data processing system,those of ordinary skill in the art will appreciate that the processes ofthe present invention are capable of being distributed in the form of acomputer readable medium of instructions, including microcode, and avariety of forms and that the present invention applies equallyregardless of the particular type of signal bearing media actually usedto carry out the distribution. Examples of computer readable mediainclude recordable-type media such a floppy disc, a hard disk drive, aRAM, and CD-ROMs and transmission-type media such as digital and analogcommunications links.

The description of the present invention has been presented for purposesof illustration and description, but is not intended to be exhaustive orlimited to the invention in the form disclosed. Many modifications andvariations will be apparent to those of ordinary skill in the art. Theembodiment was chosen and described in order to best explain theprinciples of the invention, the practical application, and to enableothers of ordinary skill in the art to understand the invention forvarious embodiments with various modifications as are suited to theparticular use contemplated.

What is claimed is:
 1. A method for controlling data transfer between amemory controller and a master device by a bus arbiter, wherein themaster device does not employ buffer reservations for data transfers,the method comprising the steps of: granting a first unrequested databus grant to a memory controller; in response to receiving a data busrequest by a master device, rescinding the first unrequested data busgrant; in response to rescinding the unrequested data bus grant,resetting a delay counter; incrementing the delay counter per cycle; andpreventing a second data bus grant while the delay is not equal to apreconfigured value.
 2. The method of claim 1 further comprising: inresponse to giving a data bus grant to the memory controller, setting apark-enabled flag.
 3. The method of claim 2 further comprising: checkingthe park-enabled flag before giving the unrequested data bus grant. 4.The method of claim 1 further comprising: in response to receiving abuffer-full signal from the master device to the bus arbiter, rescindingthe unrequested data bus grant.
 5. The method of claim 1 wherein thearbiter provides fair and unprioritized arbitration.
 6. The method ofclaim 1 wherein access to the data bus is arbitrated by an arbiter in anode controller.
 7. The method of claim 6 wherein the master device is anode controller.
 8. An apparatus for controlling data transfer between amemory controller and a master device by a bus arbiter, wherein themaster device does not employ buffer reservations for data transfers,the apparatus comprising: granting means for granting a firstunrequested data bus grant to a memory controller; rescinding means forrescinding, in response to receiving a data bus request by a masterdevice, the first unrequested data bus grant; resetting means forresetting, in response to rescinding the unrequested data bus grant, adelay counter; incrementing means for incrementing the delay counter percycle; and preventing means for preventing a second data bus grant whilethe delay is not equal to a preconfigured value.
 9. The apparatus ofclaim 8 further comprising: setting means for setting a park-enabledflag in response to giving a data bus grant to the memory controller.10. The apparatus of claim 9 further comprising: checking means forchecking the park-enabled flag before giving the unrequested data busgrant.
 11. The apparatus of claim 8 further comprising: rescinding meansfor rescinding, in response to receiving a buffer-full signal from themaster device to the bus arbiter, the unrequested data bus grant. 12.The apparatus of claim 8 wherein the arbiter provides fair andunprioritized arbitration.
 13. The apparatus of claim 8 wherein accessto the data bus is arbitrated by an arbiter in a node controller. 14.The apparatus of claim 13 wherein the master device is a nodecontroller.
 15. A computer program product in a computer-readable mediumfor use in a data processing system for controlling data transferbetween a memory controller and a master device by a bus arbiter,wherein the master device does not employ buffer reservations for datatransfers, the computer program product comprising: instructions forgranting a first unrequested data bus grant to a memory controller;instructions for rescinding, in response to receiving a data bus requestby a master device, the first unrequested data bus grant; instructionsfor resetting, in response to rescinding the unrequested data bus grant,a delay counter; instructions for incrementing the delay counter percycle; and instructions for preventing a second data bus grant while thedelay is not equal to a preconfigured value.
 16. The computer programproduct of claim 15 further comprising: instructions for setting, inresponse to giving a data bus grant to the memory controller, apark-enabled flag.
 17. The computer program product of claim 16 furthercomprising: instructions for checking the park-enabled flag beforegiving the unrequested data bus grant.
 18. The computer program productof claim 15 further comprising: instructions for rescinding, in responseto receiving a buffer-full signal from the master device to the busarbiter, the unrequested data bus grant.
 19. The computer programproduct of claim 15 wherein the arbiter provides fair and unprioritizedarbitration.
 20. The computer program product of claim 15 wherein accessto the data bus is arbitrated by an arbiter in a node controller. 21.The computer program product of claim 20 wherein the master device is anode controller.